1. Field of the Invention
This invention relates to computer programs in general, and in particular, to a method and related apparatus for maintaining coherency of secondary data in a computer system.
2. Description of the Related Art
This invention may be implemented in a wide variety of circumstances in a wide variety of computer systems. To simplify this description of the invention, however, a particular preferred embodiment of the invention will be described in detail, and comments will be added to indicate some other possible implementations or variations of the invention. Persons of skill in the art will recognize still other possible implementations and variations.
The preferred embodiment of the invention is implemented in a virtual computer system comprising multiple physical processors and multiple virtual processors, although the invention may also be advantageously implemented in a uniprocessor system. The multiple physical and virtual processors of the preferred embodiment may also or alternatively comprise multiple logical processors, such as those found in processors that include simultaneous multi-threading, which is known as “Hyper-Threading” in terminology introduced by Intel Corporation. Logical processors and simultaneous multi-threading technology will be described in greater detail below. The preferred embodiment is used to maintain TLB (Translation Lookaside Buffer) coherency in a multiprocessor virtual computer system. More specifically, the virtual computer system in which the preferred embodiment is implemented comprises one or more virtual machine monitors (VMMs) running on a physical, multiprocessor computer system, with each VMM supporting a virtual machine (VM). Accordingly, this background section provides introductions to the technologies of virtual memory systems, multiprocessor systems and simultaneous multi-threaded CPUs, TLB coherency issues and techniques, and virtual machine monitors. Next, this background section provides a description of a general virtualized computer system in which the invention may be implemented.
Virtual Memory Systems
The design and use of virtual memory systems are well known in the art, and there are numerous books and other technical references available on the subject. This invention may be implemented in various different computer systems, using various different virtual memory techniques. For purposes of an example only, the invention will be described in relation to a virtual memory system based on the x86 architecture from Intel Corporation. The invention also applies to other computer systems with other address translation mechanisms. This architecture is described in the IA-32 Intel Architecture Developer's Manual, a three-volume set, which is currently available on the Internet website of Intel Corporation, and which is hereby incorporated by reference. Volume 3 of that set, the Software Developer's Manual, is particularly informative regarding the virtual memory functions of the architecture.
FIG. 1 is a block diagram of the major functional components of a general virtual memory system in a computer. The system comprises a CPU 10, a memory management unit (MMU) 12, a TLB 14, a random access memory (RAM) 16, a plurality of page tables 18, an operating system (OS) 22, a software memory manager (S-MMU) 24, a hard disk drive 20 and a direct memory access (DMA) controller 21. Each of the functional units illustrated in FIG. 1 may be implemented by conventional components of the well-known personal computer (PC) standard architecture. The CPU 10 may also be called a processor. The RAM 16 may also be called a primary memory, while the hard drive 20 may be called a secondary memory. Also, the MMU 12 and the S-MMU 24 may be considered parts of a more general memory management unit, in which the S-MMU 24 is the software that controls the hardware MMU 12. The CPU 10 and the MMU 12 may be combined within a single integrated circuit (IC) component, or they may be separate components. Also, the TLB 14 may be contained within the same IC component as the MMU 12, or it may be a separate device.
The most basic function of the CPU 10 is to execute computer programs, including the OS 22. The computer programs are generally stored on the hard drive 20 and loaded into the RAM 16 for execution. The CPU 10 issues memory read commands to retrieve instructions of the computer programs from the RAM 16 and then executes the retrieved instructions. The execution of instructions requires a myriad of other functions too, including reading data from and writing data to the RAM 16. For example, an instruction executed by the CPU 10 may require an operation to be performed on an operand, which may be located in the RAM 16, or the instruction may require that a value be written to a stack, which may also be located within the RAM 16. All information stored in the RAM 16 may be called data, whether the data consists of instructions, operands, stack data or other types of data. At times, however, a distinction may be drawn between different types of data. In addition, the term “computer program” will generally include instructions, operands and the associated stack.
A computer program is loaded from the hard drive 20 into the RAM 16 for execution because fetching information from the RAM 16 is much quicker than from the hard drive 20, which enables the CPU 10 to execute the program much more quickly. Earlier computer systems would load an entire computer program into the RAM 16 for execution, including providing additional RAM required by the program during execution, such as for a data stack. However, RAM is relatively expensive in comparison to the cost of other data storage devices, such as disk drives. As a result, computer systems are often designed with a limited amount of RAM, in comparison to the address space of the system, especially in systems that use 64-bit addressing. This gives rise to various situations in which a computer program requires more memory space than is available in the RAM 16. A simple example of such a situation is when a computer program is simply larger than the RAM 16 of the system on which the program is to run. Another example is in a multiprocessing system, when the sum of the memory required by all of the executing threads and the OS 22 exceeds the amount of RAM 16 in the computer system. Virtual memory techniques may be used to enable the execution of a computer program in such a situation where the RAM 16 that is available for use is less than the total amount of memory required by a computer program.
Virtual memory techniques may be implemented, for example, using a combination of hardware and software. The software portion of such an implementation may be provided by the S-MMU 24 of the OS 22 of FIG. 1, while much of the hardware functionality may be provided by the MMU 12. The MMU 12 may be included, along with the CPU 10, within a single microprocessor device, such as an Intel Pentium microprocessor, or the MMU 12 may be a separate device. Virtual memory techniques give the appearance, to a computer program, that there is more RAM available than is really the case. The computer program is provided with a virtual address space, which contains all of its instructions, data and stack. The virtual address space is generally larger than the available RAM 16, but the computer program may use the entire virtual address space as if it were all contained in the RAM 16. The virtual address space may have various different types of organization, such as linear or segmented. At any given time, one or more parts of the computer program will be in the RAM 16 while one or more other parts of the computer program will not be in the RAM 16, but will be stored on the hard drive 20. If the computer program attempts to use a part of its address space that is currently not contained in the RAM 16, the S-MMU 24 will typically transfer the required part of the computer program from the hard drive 20 to the RAM 16.
To implement a virtual memory system, a computer program may be divided into a number of units called pages. For this discussion, assume a 4 kilobyte (Kbyte) page, which is one possible page size in the x86 architecture. Some of the pages of the computer program are loaded into the RAM 16, while others are not, depending on the amount of the RAM 16 that is available to the computer program. Also, the pages that are loaded into the RAM 16 may not be loaded contiguously. Typically, a particular page of the computer program on the hard drive 20 could be loaded into any available page within the RAM 16.
During execution of a computer program or process, the CPU 10 generates addresses within the virtual address space of the computer program or process, for reading data from and writing data to the RAM 16. As used herein, a process is generally an instance of a computer program. The addresses generated by the CPU 10 may be called virtual addresses or linear addresses. However, the virtual addresses cannot be directly applied to the RAM 16 in a virtual memory system to access the desired memory locations. Instead, the virtual addresses must first be translated into corresponding physical addresses within a physical address space. The physical address space comprises the addresses that are used to access specific memory locations within the RAM 16. The MMU 12 and the S-MMU 24 have primary responsibility for translating or mapping addresses from the virtual address space to the physical address space. When the CPU 10 attempts to access data from the computer program that resides on a page of the program that is not currently loaded into the RAM 16, the MMU 12 determines that the page is not resident in the RAM 16, a page fault occurs and a trap to the OS 22 ensues. The S-MMU 24 subsequently transfers the required page from the hard drive 20 into the RAM 16. After the page transfer is complete, execution of the computer program resumes at the same instruction that resulted in the page fault. This time, however, the MMU 12 will determine that the page is loaded into the RAM 16 and the memory access will be completed successfully. If there is not enough available space in the RAM 16 for loading the required page during a page fault, the S-MMU 24 typically ejects another page from the RAM 16, and the space that the ejected page was occupying is freed up for loading the new page. If the page that is being ejected has been modified in the RAM 16 since it was loaded from the hard drive 20, then it is written back to the hard drive 20 before its memory space is used for the new page.
As described in greater detail below, the MMU 12 initially uses the page tables 18, located within the RAM 16, to translate virtual addresses into physical addresses. In this process, when the MMU 12 receives a virtual address from the CPU 10 for a memory read or write, the MMU 12 must first perform at least one memory read within the page tables 18 just to determine the corresponding physical address. The MMU 12 must then perform another memory access to complete the read or write required by the CPU 10. If the MMU 12 had to access the page tables 18 for every memory access from the CPU 10, using the virtual memory system would add at least one extra memory cycle to each memory access. In some virtual memory systems, multiple memory accesses are required to map a virtual address to a physical address, using the page tables 18. The added memory cycles would slow down the execution of instructions, which would reduce the overall processing power of the computer system. The primary purpose of the TLB 14 is to reduce the number of additional memory accesses that are required to implement the virtual memory system. The TLB 14 is basically a cache for page table entries and typically is located within the MMU 12. Fortunately, when a CPU 10 is executing a computer program, most of its memory accesses will be to a limited number of pages within the RAM 16. At any given time, for a particular program, the CPU 10 will typically access one or a few pages of code, one or a few pages of data and one or a few pages for the stack, depending on the page size used.
At this point, it is useful to discuss page numbers. As described above, the virtual address space of a computer program or a process is divided into a number of pages. Each of these pages can be numbered consecutively, resulting in virtual page numbers. In the same way, the physical address space of the RAM 16 can be divided into pages as well. These pages can also be numbered consecutively, resulting in physical page numbers. Now, a virtual address can be viewed as specifying a virtual page number in the upper bits and an offset within that page in the lower bits. In the same way, a physical address can be viewed as a physical page number combined with an offset into that physical page. For example, in a system having 32-bit addresses and a 4 Kbyte page size, such as an x86 system, the upper 20 bits of an address can be viewed as a page number and the lower 12 bits can be viewed as an offset within a given page. Then, so long as both virtual pages and physical pages begin at an address that is a multiple of the 4 Kbyte page size, the address translation process can be viewed as converting the upper address bits from a virtual page number to a physical page number, with the lower address bits remaining unchanged as the offset into the respective pages.
The MMU 12 uses the page tables 18 to perform this translation from virtual page numbers to physical page numbers. When the MMU 12 receives a virtual address from the CPU 10, the MMU 12 reads the virtual page number from the upper address bits of the address. The MMU 12 then reads information from the page tables 18 relating to the desired virtual page number. First, the page tables 18 will indicate whether the virtual page number is currently loaded into the RAM 16. If the virtual page is not loaded into the RAM 16, a page fault is generated and the required virtual page is loaded into the RAM 16 as described above. If the virtual page is loaded into the RAM 16, the page tables 18 will also indicate the physical page number that corresponds to the virtual page number. The MMU 12 then uses the retrieved physical page number, along with the offset from the virtual address to access the desired location within the RAM 16. In addition, the MMU 12 writes the virtual page number and the physical page number into an entry in the TLB 14, indicating the mapping between the pages. In other systems, the S-MMU 24 may be responsible for loading such mappings into the TLB 14. Accessing the page tables 18 in this manner to determine a mapping from a virtual page number to a physical page number is called walking the page tables 18. Now that the mapping from the virtual page number to the physical page number has been written into the TLB 14, if a subsequent memory access is to the same virtual page number, the MMU 12 can find the appropriate mapping in the TLB 14 within the MMU 12, without having to access the page tables 18 in the RAM 16.
The MMU 12 is designed such that the access to the TLB 14 is much quicker than an access to the page tables 18. The TLB 14 can typically only hold a relatively small number of page mappings, such as 8 to 64 entries, in comparison to the size of the page tables 18. As a result, entries must be evicted from the TLB 14 from time to time. Typically, when the MMU 12 walks the page tables 18 to determine a new mapping, the MMU 12 will evict an existing entry in the TLB 14 to make space to enter the new mapping. Thus, when the MMU 12 receives a virtual address from the CPU 10, the MMU 12 may first access the TLB 14 to determine if the desired mapping is there. If the mapping is not in the TLB 14, then the MMU 12 must perform a page table walk, as described above and in greater detail below.
FIG. 2A shows a 32-bit virtual address 30, comprising a 10-bit page directory entry (PDE) 30A, a 10-bit page table entry (PTE) 30B and a 12-bit offset 30C. FIG. 2B illustrates the structure and operation of the page tables of the x86 architecture, as a more detailed example. FIG. 2B also shows a page directory 40 with 1024 page table base address (PTBA) entries 42, including one specific PTBA entry 42X. FIG. 2B also shows a plurality of page tables 50A, 50X and 50N. These page tables, along with other page tables that are not illustrated, will be collectively referred to as page tables 50. This convention, of using a common numeric portion to refer collectively to all items having alphanumeric references containing the same numeric portion, is used throughout this description. As shown relative to the page table 50X, each of the page tables 50 comprises 1024 physical page base address (PPBA) entries 52. Page table 50X includes one specific PPBA entry 52X. FIG. 2B also shows a plurality of physical pages 60, including the physical pages 60A, 60X and 60N. As shown relative to the physical page 60X, each of the physical pages 60 comprises 4096 addressable bytes 62. Physical page 60X includes one specific byte 62X. Each of the physical pages 60, the page tables 50 and the single page directory 40 reside in the RAM 16. Each of the physical pages 60 includes 4096 bytes, or 4 Kbytes. As described above, the physical pages 60 and the virtual pages of the example in this description include 4 Kbytes of data. Each of the 1024 PTBA entries 42 in the page directory 40 comprises 32 bits, or 4 bytes. Thus, the page directory 40 also constitutes a full 4 Kbyte page in the RAM 16. Each of the 1024 PPBA entries 52 in the page tables 50 also comprises 32 bits. So, each of the page tables 50 also constitutes a full 4 Kbyte page in the RAM 16.
When the MMU 12 receives a virtual address 30 from the CPU 10, the MMU 12 may first check to see if there is an entry in the TLB 14 that provides a mapping from the virtual page number to a corresponding physical page number. The combination of the PDE 30A and the PTE 30B is considered the virtual page number 30AB. In this architecture, the TLB 14 maps 20-bit virtual page numbers to 20-bit physical page numbers. So, the MMU 12 checks whether there is a valid entry in the TLB 14 matching the virtual page number 30AB. If there is, the MMU 12 uses this entry to obtain the desired mapping to a physical page 60. If there is no matching entry in the TLB 14, the MMU 12 must walk the page tables 18. In the x86 architecture, the page directory 40 may be considered a page table 18, as well as the page tables 50. To walk the page tables 18, the MMU 12 first reads a 20-bit value from a control register CR3 in some modes of operation. This 20-bit value is used as the upper 20 bits of a 32-bit address that points to the base of the page directory 40. The lower 12 bits of this address are set to zero. Thus, the page directory 40 must begin at an address that is a multiple of the 4 Kbyte page size. The page tables 50 and the physical pages 60 must also begin at an address that is a multiple of the 4 Kbyte page size for the same reason.
Once the base address of the page directory 40 is determined, the PDE 30A is used as an index into the 1024-entry table of the page directory 40. More specifically, the 20 bits from the control register CR3 are used as the upper address bits, the 10 bits from the PDE 30A are used as the next lower address bits, and the last two address bits are set to 0 to form a memory address, which addresses the PTBA entry 42X. As illustrated in FIG. 2B, the control register CR3 points to the beginning of the page directory 40, while the PDE 30A points to the PTBA entry 42X. One bit of the PTBA entry 42X indicates whether the PTBA entry 42X is a valid entry. If it is not a valid entry, a page fault results, which may indicate an error condition in the S-MMU 24 or in the running application. If the entry is valid, a 20-bit value from the PTBA entry 42X is used as the upper bits of a base address for the page table 50X. The PTE 30B is used as an index into the 1024-entry table of the page table 50X. As shown in FIG. 2B, the page table base address entry 42X points to the base of the page table 50X, while the PTE 30B points to the PPBA entry 52X. One bit of the PPBA entry 52X indicates whether the virtual page number 30AB is currently loaded into the RAM 16. If the virtual page number 30AB is not currently loaded into the RAM 16, a page fault results and the required virtual page is loaded into the RAM 16, as described above. If the virtual page number 30AB is loaded into the RAM 16, a 20-bit value from the PPBA entry 52X is used as the upper address bits of a base address for the page table 60X for the current memory access. The offset 30C is now used as an index into the physical page 60X to identify a specific byte address 62X for the memory access. In other words, the 20 bits from the PPBA entry 52X are combined with the 12 bits from the offset 30C to form a 32-bit physical address that is used to perform the memory access requested by the CPU 10. As shown in FIG. 2B, the PPBA entry 52X points to the base of the physical page 60X, while the offset 30C points to the required byte address 62X for the memory access.
Generally, the S-MMU 24 of the OS 22 is responsible for creating and maintaining the page tables 18 for the use of the MMU 12. Either the MMU 12 or the S-MMU 24 is generally responsible for loading values into the TLB 14 for recently obtained mappings between virtual page numbers and physical page numbers. Values may be flushed from the TLB 14 either by the MMU 12 or by the S-MMU 24, or possibly by other software within the RAM 16, such as user-level application programs. Each entry within a page table 18 generally contains, in addition to a physical page number, a few other bits for indicating whether the entry is valid, what types of access are allowed for the page, whether the page has been modified and/or referenced since it was loaded into the RAM 16 and whether caching is disabled for the page. An entry within the TLB 14 generally contains a virtual page number and a physical page number, as well as a few additional bits to indicate whether the entry is valid, whether the page has been modified since being loaded into the RAM 16 and what types of access are allowed for the page. When a memory access is performed, if the MMU 12 determines that the virtual page is loaded into the RAM 16, the MMU 12 also accesses these additional bits of the entry within either the page tables 18 or the TLB 14, to determine if the requested memory access is permitted. For example, the access bits may indicate that only read accesses are permitted. If the CPU 10 attempts to write data to such a location, the MMU 12 will generate a page fault.
When a mapping for a particular virtual page number is not contained within the TLB 14 and a page table walk is performed, the MMU 12 typically evicts an entry from the TLB 14 to free up space for a new entry for the current mapping. The virtual page number will be written into the newly available entry in the TLB 14, along with the physical page number that was just determined. The additional bits within the entry of the TLB 14 are typically copied from the corresponding additional bits in the corresponding page table entry. When an entry in the TLB 14 is evicted, a bit indicating whether the page has been modified is typically copied from the entry of the TLB 14 to the corresponding entry in the page table 18. Also, if the S-MMU 24 removes a virtual page from the RAM 16 for which there is an entry in the TLB 14, the S-MMU 24 must modify the entry in the TLB 14 to indicate that the mapping is no longer valid. Other programs may also be allowed to indicate that an entry of the TLB 14 is invalid, including possibly user-level applications. The x86 architecture provides an instruction, Inylpg(virtual address), for this purpose, although the Inylpg instruction is usually restricted for use only by kernel level code. The x86 architecture is defined such that, if an entry in the TLB 14 is set as invalid, the MMU 12 will walk the page tables to determine a mapping for the virtual address. However, if an entry in the TLB 14 is not set as invalid, the MMU 12 may use the TLB 14 to obtain a mapping, or the MMU 12 may walk the page tables to determine the mapping. The x86 architecture also provides an instruction for flushing the entire contents of the TLB 14. As described above, entries within the TLB 14 may also be evicted by the MMU 12 to free up space for a new mapping for a new virtual address. Thus, an entry in the TLB 14 may be created for a specific virtual page number in response to a first access to that virtual page. During a subsequent access to the same virtual page, if the entry in the TLB 14 has been evicted by the MMU 12 in between the two accesses, a page table walk will nonetheless be required. This situation is described as a leakage of the TLB 14.
Multiprocessor Systems and Simultaneous Multi-Threaded CPUs
Multiprocessor systems are also well known in the art and there are various architectures currently available. FIG. 3 illustrates one general architecture for a multiprocessor system. FIG. 3 shows a shared primary memory 16B, an OS 22B, an S-MMU 24B, a plurality of page tables 18B, and a shared secondary memory 20B. These functional units perform the same basic functions as the corresponding functional units shown in FIG. 1, but they may need to be modified to perform these functions in a multiprocessor environment. There are various types of operating systems 22B for use in multiprocessor systems. Some multiprocessor operating systems 22B use a single operating system image to manage the entire set of processors in concert. In other multiprocessor systems, the system hardware provides a physical partitioning of the system, allowing a different instance of a multiprocessor operating system 22B to manage each partition. In the case of a multiprocessor OS 22B comprising a separate instance for each CPU, the separate instances of the OS 22B may be executed in separate private memories associated with each of the multiple CPUs. The shared primary memory 16B may be the same type of memory as the RAM 16, except perhaps larger, and the shared secondary memory 20B may be the same as the hard drive 20, except perhaps larger. The page tables 18B may be the same as the page tables 18, except that there may be more sets of page tables 18B because of the multiple CPUs.
FIG. 3 also shows a first processor (such as a microprocessor) 9A, having a first CPU 11A, a first MMU 13A and a first TLB 15A. The microprocessor 9A is also connected to a first private memory 17A. FIG. 3 also shows a second processor (such as a microprocessor) 9B, having a second CPU 11B, a second MMU 13B and a second TLB 15B. The microprocessor 9B is also connected to a second private memory 17B. The multiprocessor system of FIG. 3 may also have additional microprocessors 9 and associated private memories 17. The microprocessors 9 may be, for example, based on the x86 architecture. In addition, each of the private memories 17 is optional.
In a single-processor system, there may be a single set of page tables 18 or there may be multiple sets of page tables 18. Each process could have its own set of page tables 18 or there could be some sharing of page tables 18. In a multiprocessor system, there could be page tables 18B in the shared primary memory 16B, in one or more of the private memories 17, or both, and any of these page tables 18B could be shared between multiple processes or exclusive to a single process. As another alternative to the system illustrated in FIG. 3, one or more TLBs 15 could be shared among multiple CPUs 11. One example of a multiprocessor system having a shared TLB 15 is illustrated in FIG. 4 and described below.
The virtual memory system implemented in the system of FIG. 3 can be functionally similar to the virtual memory system described above in connection with FIGS. 1 and 2. More specifically, the TLBs 15 and the page tables 18B can have the same basic structure and functionality as the TLB 14 and the page tables 18, respectively, and the MMUs 13 and the S-MMU 24B can control and use the TLBs 15 and the page tables 18B in the same general manner that the MMU 12 and the S-MMU 24 control and use the TLB 14 and the page tables 18. If there is no sharing of the TLBs 15 or the page tables 18B between multiple CPUs 11, then the virtual memory system of FIG. 3 can be functionally the same as the virtual memory system of FIGS. 1 and 2, but with a separate instance of the virtual memory system for each of the CPUs 11. However, if there is any sharing of the TLBs 15 or the page tables 18B between the multiple CPUs 11, the virtual memory system gets more complicated. The following discussion will focus on a multiprocessor system containing only two CPUs 11, for simplicity, although it also applies to systems with more CPUs 11.
The discussion also applies to systems that have only one physical CPU, if the CPU implements simultaneous multi-threading techniques. Simultaneous multi-threading techniques are known in the art and are becoming more prevalent, especially in high-performance CPUs, such as the Xeon microprocessor from Intel Corporation. In a CPU that implements simultaneous multi-threading, multiple instruction streams are executed simultaneously. With multiprogramming or multithreading techniques, in contrast, different instruction streams are executed during separate time slices. A CPU that does not provide simultaneous multi-threading can generally be modeled as an interpreter loop, in which the CPU repeatedly fetches an instruction, fetches any required operands, performs an operation and does something with the result of the operation, such as writing the result to memory, before moving on to fetch the next instruction. A simultaneous multi-threaded CPU, in contrast, can be modeled as multiple independent interpreter loops running concurrently. Effectively, the single physical CPU core provides the capabilities of multiple logical CPUs. However, a simultaneous multi-threaded processor may have only a single TLB, or it may have multiple TLBs. For the purposes of this invention and the discussion below, a simultaneous multi-threaded processor, having multiple logical CPUs but only one TLB, is functionally equivalent to a multiprocessor system having multiple physical CPUs and a single, shared TLB. This invention and the following discussion may apply to any computer system in which multiple processes or threads are executing simultaneously on multiple physical or logical CPUs, and the multiple processes or threads share a common TLB or page table. In fact, as will become apparent below, this invention may even apply in a system having a separate TLB and a separate set of page tables for each process or thread, so long as at least one process or thread has write access to the page tables of at least one other process or thread, even if such access is provided inadvertently, such as due to a system software error, for example.
FIG. 4 illustrates another example architecture for a multiprocessor computer system. Specifically, FIG. 4 shows a first CPU 11C, a second CPU 11D, an MMU 13C, a shared TLB 15C, a shared primary memory 16C, an OS 22C, an S-MMU 24C, a set of page tables 18C, and a shared secondary memory 20C. Each of the CPUs 11C and 11D may be either physical or logical, and there may also be additional physical and/or logical CPUs 11. The MMU 13C and the TLB 15C are shared between the CPUs 11C and 11D. Otherwise, the functional units illustrated in FIG. 4 may be equivalent to the corresponding functional units illustrated in FIG. 3.
TLB Coherency Issues and Techniques
Referring again to the multiprocessor system of FIG. 3, suppose that the CPU 11A is executing a first process and the CPU 11B is executing a second process. The system of FIG. 3 implements a virtual memory system, with some pages of the virtual address space of the first process loaded into primary memory 16B and others remaining in the secondary memory 20B. Suppose, for the moment, that the first and second processes share a common set of page tables 18B. The page tables 18B indicate, for each virtual page, whether it is loaded into the primary memory 16B or whether it remains in the secondary memory 20B. The page tables 18B also indicate, for each virtual page loaded into the primary memory 16B, the corresponding physical page number into which the virtual page is loaded. The TLB 15A may also contain one or more entries indicating mappings between virtual pages and physical pages of the primary memory 16B.
Suppose further that the first process executes a first instruction that accesses a first memory location on a first virtual page that is currently loaded into a first physical page. Suppose that the MMU 13A walks the page tables 18B to determine a mapping between the first virtual page and the first physical page, and stores this mapping in the TLB 15A. Now suppose that the second process changes the page tables 18B, or performs some action that causes the page tables 18B to be changed. For example, the second process may attempt to access a second virtual page that is not currently loaded into the primary memory 16B, causing a page fault. In response to the page fault, the S-MMU 24B loads the second virtual page from the secondary memory 20B into the primary memory 16B. Suppose that the S-MMU 24B loads the second virtual page from the secondary memory 20B into the first physical page of the primary memory 16B, replacing the first virtual page. The S-MMU 24B updates the page tables 18B to indicate that the second virtual page is now loaded into the primary memory 16B and is mapped to the first physical page, and to indicate that the first virtual page is no longer loaded into the primary memory 16B.
Now suppose that the first process executes a second instruction that again accesses the first memory location on the first virtual page, or some other memory location on the first virtual page. If the MMU 13A accesses the TLB 15A to determine a mapping for the first virtual page, the previously stored mapping will indicate that the first virtual page is mapped to the first physical page. The MMU 13A would then retrieve the contents of the corresponding memory location within the first physical page and provide this data to the CPU 11A for executing the second instruction. However, the data retrieved by the MMU 13A is actually from the first physical page, instead of from the second physical page as indicated in the page tables. Thus, the CPU 11A would execute the second instruction based on incorrect data, possibly corrupting the data for the first process, the second process, or both. If the first virtual page contained code for the first process, as opposed to operand data or stack data, so that the attempted memory access were an instruction fetch, then the CPU 11A would attempt to execute whatever data is retrieved by the MMU 13A. If the second virtual page happens to contain operand or stack data, then the CPU 11A would nonetheless attempt to interpret the returned data as an instruction and try to execute the interpreted instruction. This situation would also likely lead to corrupted data, or worse.
This situation is referred to as an incoherent TLB. The TLB 15A operates as a cache for the virtual addressing data contained in the page tables 18B. If an entry in a page table 18B is loaded into the TLB 15A, and then the same entry in the page table 18B is modified, then the corresponding entry in the TLB 15A should generally be updated accordingly to maintain coherency with the page table 18B, or at least invalidated to avoid an incoherency.
Multiprocessor systems, as well as uniprocessor systems, generally provide methods to avoid TLB incoherencies. One common technique in a multiprocessor system, referred to as a “TLB shootdown,” enables the OS 22B to cause the mapping between the first virtual page and the first physical page in the TLB 15A to be flushed or invalidated. First, the OS 22B, executing on the CPU 11B, may coordinate with the CPU 11A to bring the CPU 11A into a safe state, so that the CPU 11A is not using the first virtual page or the page table entry mapping the first virtual page to the first physical page, although this step may be omitted in many cases. Next, the OS 22B, again executing on the CPU 11B, switches the first physical page to contain the second virtual page instead of the first virtual page and modifies the appropriate entries in the page table 18B. Next, the OS 22B causes the CPU 11A to flush or invalidate the entry in the TLB 15A that maps the first virtual page to the first physical page. In the x86 architecture, for example, the OS 22B may cause the CPU 11A to execute an Inylpg instruction, specifying the page number for the first virtual page. Now, when the second instruction is executed by the CPU 11A, causing the second access to the first memory location, the MMU 13A does not find a mapping for the first virtual page in the TLB 15A and is forced to walk the page tables 18B. The MMU 13A now determines that the first virtual page is no longer loaded into the primary memory 16B, as appropriate.
Virtual Machine Monitors
A virtual machine monitor (VMM) is a piece of software that runs on top of the hardware of a computer system having a first hardware platform and creates an abstracted or virtualized computer system having a second hardware platform. The second hardware platform, or virtualized platform, may be the same as, similar to, or substantially different from, the first hardware platform. The VMM exports all of the features of the virtualized platform, to create a virtual machine (VM) that is functionally equivalent to an actual hardware system implementing the second hardware platform. The VMM generally performs all of the functions that would be performed by a physical implementation of the virtualized hardware platform, to achieve the same results. For example, a VMM generally implements a virtual memory system that is functionally equivalent to the virtual memory system that would result from a physical implementation of the virtualized platform. Various designs for such VMMs are well known in the art, and this invention may be implemented in any such design.
An OS designed to run on a computer system having the virtualized hardware platform can be loaded on top of the VMM, and the OS should not be able to determine that it is not running directly on an actual hardware system implementing the virtualized hardware platform. Therefore, in the case where the virtualized hardware platform is the same as the physical hardware platform, the OS can be loaded directly onto the actual computer system or on top of the VMM, and the OS would not be able to determine whether the machine on which it is running is the physical machine or the virtual machine. Drivers and other system software that are designed for computer systems having the virtualized hardware platform can also be loaded onto the VMM. An OS running on a VMM, along with drivers and other system software, is called a guest OS. In addition, application programs that are designed to operate on the guest OS may also be loaded onto the VMM. An application program loaded onto the VMM is called a guest application.
As one example of a VMM implementation, a VMM may run on an x86 computer system, and it may virtualize an x86 system. In this case, the VMM creates a VM that is compatible with the x86 architecture. Any operating system that can run on an x86 system may be loaded on top of the VMM. For example, a Windows OS from Microsoft Corporation, such as the Windows XP OS, may be loaded as the guest OS on top of the VMM. Application programs that are designed to operate on a system running the Windows XP OS can then also be loaded onto the VMM. The guest OS and the application programs will execute just as if they were loaded directly onto the underlying physical x86 system.
VMMs have also been designed to operate on multiprocessor systems, and to virtualize multiprocessor systems. For example, one or more VMMs may execute on the hardware platform illustrated in FIG. 3, and may create a VM having the same hardware architecture. As described above, the VMMs should generally be functionally equivalent to the virtualized hardware platform. Of particular relevance to this invention, the VMMs should generally provide a virtual memory system that is functionally equivalent to the virtual memory system of the virtualized platform. Thus, the VMMs should virtualize one or more MMUs and one or more TLBs that are functionally equivalent to the MMUs and TLBs of an actual physical implementation of the virtualized hardware platform. The virtualized MMUs and TLBs should interact with the S-MMU of the guest OS and with the guest applications in the same manner as the corresponding physical MMUs and TLBs would interact with the S-MMU and the guest applications. In particular, the VMMs should provide the same safeguards against virtual memory conflicts, which could result from shared TLBs, shared page tables and/or a shared primary memory, as are provided by a physical implementation of the virtualized platform. For example, suppose a VMM, or a set of VMMs, were designed to virtualize the hardware platform illustrated in FIG. 3, including the function described above in which the second CPU 11B is able to initiate a TLB shootdown to cause the first CPU 11A to flush or invalidate a mapping in the first TLB 15A. Such a VMM would virtualize the first microprocessor 9A, the second microprocessor 9B and the first TLB 15A such that the second virtual CPU could communicate with the first virtual CPU causing the first virtual CPU to flush the mapping in the virtual TLB.
General Virtualized Computer System
FIG. 5 illustrates, in part, the general configuration of a virtual computer system 700, including a virtual machine 200, which is installed as a “guest” on a “host” hardware platform 100. As FIG. 5 shows, the hardware platform 100 includes one or more processors (CPUs) 110, system memory 130, and one or more local storage devices, which typically includes a local disk 140. The system memory is typically some form of high-speed RAM, whereas the disk (one or more) is typically a non-volatile, mass storage device. The hardware 100 also typically includes other conventional mechanisms such as a memory management unit (MMU) 150 and various registers 160. The hardware 100 may also include one or more interface cards for interfacing with external devices, computers, systems and/or networks. For example, the hardware 100 may include a data interface 170 for connecting to an external data storage system or network and/or a network interface 180 for connecting to a computer network.
Each VM 200 typically includes at least one virtual CPU 210, at least one virtual disk 240, a virtual system memory 230, a guest operating system 220 (which may simply be a copy of a conventional operating system), and various virtual devices 280, in which case the guest operating system (“guest OS”) includes corresponding drivers 224. All of the components of the VM may be implemented in software using known techniques to emulate the corresponding components of an actual computer.
One or more guest applications 260 may also be loaded into the VM 200. Executable files will be accessed by the guest OS from a virtual disk or virtual memory, which may simply be portions of an actual physical disk or memory allocated to that VM. Once an application is installed within the VM, the guest OS retrieves files from the virtual disk just as if they had been pre-stored as the result of a conventional installation of the application. The design and operation of virtual machines is well known in the field of computer science.
A VMM, or some such interface, is generally required between a VM and the underlying host platform (in particular, the CPU), which is responsible for actually executing VM-issued instructions and transferring data to and from the actual memory and storage devices. In FIG. 5, a VMM is shown as component 300. A VMM is usually a thin piece of software that runs directly on top of a host, or directly on the hardware, and virtualizes the resources of the physical host machine. Among other components, the VMM therefore usually includes device emulators 330, which may constitute the virtual devices 280 that the VM 200 accesses.
The VMM also usually tracks and either forwards (to some form of operating system) or itself schedules and handles all requests by its VM for machine resources, as well as various faults and interrupts. A mechanism known in the art as an exception or interrupt handler 355 is therefore included in the VMM. As is well known, such an interrupt/exception handler normally includes an interrupt descriptor table (IDT), or some similar table, which is typically a data structure that uses information in the interrupt signal to point to an entry address for a set of instructions that are to be executed when the interrupt/exception occurs. FIG. 5 also shows, within the VMM 300, a patch producer 396 and a patch consumer 399, which will be described below in connection with the preferred embodiment of the invention.
Although the VM (and thus the user of applications running in the VM) cannot usually detect the presence of the VMM, the VMM and the VM may be viewed as together forming a single virtual computer. They are shown in FIG. 5 as separate components for the sake of clarity. Moreover, the various virtualized hardware components such as the virtual CPU(s) 210, the virtual memory 230, the virtual disk 240, and the virtual device(s) 280 are shown as being part of the VM 200 for the sake of conceptual simplicity—in actual implementations these “components” are usually constructs or emulations exported to the VM by the VMM. For example, the virtual disk 240 is shown as being within the VM 200. This virtual component, which could alternatively be included among the virtual devices 280, may in fact be implemented as one of the device emulators 330 in the VMM. The device emulators 330 emulate the system resources for use within the VM. These device emulators will then typically also handle any necessary conversions between the resources as exported to the VM and the actual physical resources.
As in most modern computers and as described above, the address space of the memory 130 is partitioned into pages (for example, in the Intel x86 architecture) or other analogous units. Applications then address the memory 130 using virtual addresses (VAs), which include virtual page numbers (VPNs). The VAs are then mapped to physical addresses (PAs) that are used to address the physical memory 130. VAs and PAs have a common offset from a base address, so that only the VPN needs to be converted into a corresponding physical page number (PPN). Similar mappings are used in other architectures where relocatability is possible.
An extra level of addressing indirection is typically implemented in virtualized systems in that a VPN issued by an application 260 in the VM 200 is remapped twice in order to determine which page of the hardware memory is intended. The first mapping is provided by a mapping module within the guest OS 220, which translates the guest VPN (GVPN) into a corresponding guest PPN (GPPN) in the conventional manner. The guest OS therefore “believes” that it is directly addressing the actual hardware memory, but in fact it is not.
Of course, a valid address to the actual hardware memory must ultimately be generated. A memory management module 350, located typically in the VMM 300, therefore performs the second mapping by taking the GPPN issued by the guest OS 220 and mapping it to a hardware (or “machine”) page number PPN that can be used to address the hardware memory 130. This GPPN-to-PPN mapping may instead be done in the main system-level software layer (such as in a mapping module in a kernel 600, which is described below), depending on the implementation. From the perspective of the guest OS, the GVPN and GPPN might be virtual and physical page numbers just as they would be if the guest OS were the only OS in the system. From the perspective of the system software, however, the GPPN is a page number that is then mapped into the physical memory space of the hardware memory as a PPN.
In some systems, such as the Workstation product of VMware, Inc., of Palo Alto, Calif., the VMM is co-resident at system level with a host operating system. Both the VMM and the host OS can independently modify the state of the host processor, but the VMM calls into the host OS via a driver and a dedicated user-level application to have the host OS perform certain I/O operations on behalf of the VM. The virtual computer in this configuration is thus fully hosted in that it runs on an existing host hardware platform and together with an existing host OS.
In other implementations, a dedicated kernel takes the place of and performs the conventional functions of the host OS, and virtual computers run on the kernel. FIG. 5 illustrates a kernel 600 that serves as the system software for several VM/VMM pairs 200/300, . . . , 200N/300N. Compared with a system in which VMMs run directly on the hardware platform, use of a kernel offers greater modularity and facilitates provisioning of services that extend across multiple VMs (for example, for resource management). Compared with the hosted deployment, a kernel may offer greater performance because it can be co-developed with the VMM and be optimized for the characteristics of a workload consisting of VMMs. The ESX Server product of VMware, Inc., has such a configuration.
A kernel-based virtualization system of the type illustrated in FIG. 5 is described in U.S. patent application Ser. No. 09/877,378 (“Computer Configuration for Resource Management in Systems Including a Virtual Machine”), now U.S. Pat. No. 6,961,941, which is incorporated here by reference. The main components of this system and aspects of their interaction are, however, outlined below.
At boot-up time, an existing operating system 420 may be at system level and the kernel 600 may not yet even be operational within the system. In such case, one of the functions of the OS 420 may be to make it possible to load the kernel 600, after which the kernel runs on the native hardware 100 and manages system resources. In effect, the kernel, once loaded, displaces the OS 420. Thus, the kernel 600 may be viewed either as displacing the OS 420 from the system level and taking this place itself, or as residing at a “sub-system level.” When interposed between the OS 420 and the hardware 100, the kernel 600 essentially turns the OS 420 into an “application,” which has access to system resources only when allowed by the kernel 600. The kernel then schedules the OS 420 as if it were any other component that needs to use system resources.
The OS 420 may also be included to allow applications unrelated to virtualization to run; for example, a system administrator may need such applications to monitor the hardware 100 or to perform other administrative routines. The OS 420 may thus be viewed as a “console” OS (COS). In such implementations, the kernel 600 preferably also includes a remote procedure call (RPC) mechanism to enable communication between, for example, the VMM 300 and any applications 430 installed to run on the COS 420.
In kernel-based systems such as the one illustrated in FIG. 5, there must be some way for the kernel 600 to communicate with the VMM 300. In general, the VMM 300 can call into the kernel 600 but the kernel cannot call directly into the VMM. The conventional technique for overcoming this is for the kernel to post “actions” (requests for the VMM to do something) on an action queue stored in memory 130. As part of the VMM code, the VMM looks at this queue periodically, and always after it returns from a kernel call and also before it resumes a VM. One typical action is the “raise interrupt” action: If the VMM sees this action it will raise an interrupt to the VM 200 in the conventional manner.
As is known, for example, from U.S. Pat. No. 6,397,242 (Devine, et al., 28 May 2002), some virtualization systems allow VM instructions to run directly (in “direct execution”) on the hardware CPU(s) when possible. When necessary, however, VM execution is switched to the technique known as “binary translation,” during which the VM is running in the VMM. In any systems where the VM is running in direct execution when it becomes necessary for the VMM to check actions, the kernel must interrupt the VMM so that it will stop executing VM instructions and check its action queue. This may be done using known programming techniques.
The kernel 600 handles not only the various VMM/VMs, but also any other applications running on the kernel, as well as the COS 420, as entities that can be separately scheduled on the hardware CPU(s) 110. In this disclosure, each schedulable entity is referred to as a “world,” which contains a thread of control, an address space, machine memory, and handles to the various device objects that it is accessing. Worlds are stored in a portion of the memory space controlled by the kernel. More specifically, the worlds are controlled by a world manager, represented in FIG. 5 within the kernel 600 as module 612. Each world also has its own task structure, and usually also a data structure for storing the hardware state currently associated with the respective world.
The kernel 600 includes a memory management unit 616 that manages all machine memory that is not allocated exclusively to the COS 420. When the kernel 600 is loaded, the information about the maximum amount of memory available on the machine is available to the kernel, as well as information about how much of it is being used by the COS. Part of the machine memory is used for the kernel 600 itself and the rest is used for the virtual machine worlds.
Virtual machine worlds use machine memory for two purposes. First, memory is used to back portions of each world's memory region, that is, to store code, data, stacks, etc., in the VMM page table. For example, the code and data for the VMM 300 is backed by machine memory allocated by the kernel 600. Second, memory is used for the guest memory of the virtual machine. The memory management module may include any algorithms for dynamically allocating memory among the different VM's 200.
Interrupt and exception handling is related to the concept of “worlds” described above. As mentioned above, one aspect of switching worlds is changing various descriptor tables. One of the descriptor tables that is loaded when a new world is to be run is the new world's IDT. The kernel 600 therefore preferably also includes an interrupt/exception handler 655 that is able to intercept and handle (using a corresponding IDT in the conventional manner) interrupts and exceptions for all devices on the machine. When the VMM world is running, whichever IDT was previously loaded is replaced by the VMM's IDT, such that the VMM will handle all interrupts and exceptions.
The VMM will handle some interrupts and exceptions completely on its own. For other interrupts/exceptions, it will be either necessary or at least more efficient for the VMM to call the kernel to have the kernel either handle the interrupts/exceptions itself, or to forward them to some other sub-system such as the COS. One example of an interrupt that the VMM can handle completely on its own, with no call to the kernel, is a check-action IPI (inter-processor interrupt). One example of when the VMM preferably calls the kernel, which then forwards an interrupt to the COS, would be where the interrupt involves devices such as a mouse, which is typically controlled by the COS. The VMM may forward still other interrupts to the VM.
In the preferred embodiment of the invention, the kernel 600 is responsible for providing access to all devices on the physical machine. In addition to other modules that the designer may choose to load onto the system for access by the kernel, the kernel will therefore typically load conventional drivers as needed to control access to devices. Accordingly, FIG. 5 shows a module 610 containing loadable kernel modules and drivers. The kernel 600 may interface with the loadable modules and drivers in a conventional manner, using an application program interface (API) or similar interface.